L2 cache controller with slice directory and unified cache structure

ABSTRACT

A cache memory logically partitions a cache array having a single access/command port into at least two slices, and uses a first cache directory to access the first cache array slice while using a second cache directory to access the second cache array slice, but accesses from the cache directories are managed using a single cache arbiter which controls the single access/command port. In the illustrative embodiment, each cache directory has its own directory arbiter to handle conflicting internal requests, and the directory arbiters communicate with the cache arbiter. An address tag associated with a load request is transmitted from the processor core with a designated bit that associates the address tag with only one of the cache array slices whose corresponding directory determines whether the address tag matches a currently valid cache entry. The cache array may be arranged with rows and columns of cache sectors wherein a given cache line is spread across sectors in different rows and columns, with at least one portion of the given cache line being located in a first column having a first latency and another portion of the given cache line being located in a second column having a second latency greater than the first latency. The cache array outputs different sectors of the given cache line in successive clock cycles based on the latency of a given sector.

CROSS-REFERENCE TO RELATED APPLICATIONS

This application is related to U.S. patent application Ser. No.11/054,930, now U.S. Pat. No. 7,366,841 entitled “L2 CACHE ARRAYTOPOLOGY FOR LARGE CACHE WITH DIFFERENT LATENCY DOMAINS” filedconcurrently herewith, U.S. patent application Ser. No. 11/054,925entitled “SYSTEM BUS STRUCTURE FOR LARGE L2 CACHE ARRAY TOPOLOGY WITHDIFFERENT LATENCY DOMAINS” filed concurrently herewith, and U.S. patentapplication Ser. No. 11/055,262, now U.S. Pat. No. 7,308,537 entitled“HALF-GOOD MODE FOR LARGE L2 CACHE ARRAY TOPOLOGY WITH DIFFERENT LATENCYDOMAINS” filed concurrently herewith, each of which is herebyincorporated.

BACKGROUND OF THE INVENTION

1. Field of the Invention

The present invention generally relates to computer systems, and moreparticularly to a memory hierarchy for a computer system that includeslarge cache structures having different latencies across the cachearrays.

2. Description of the Related Art

The basic structure of a conventional computer system includes one ormore processing units which are connected to various peripheral devices(including input/output devices such as a display monitor, keyboard, andpermanent storage device), a memory device such as random access memory(RAM) that is used by the processing units to carry out programinstructions and store operand data, and firmware which seeks out andloads an operating system from one of the peripherals (usually thepermanent memory device) whenever the computer is first turned on. Theprocessing units typically communicate with the peripheral devices bymeans of a generalized interconnect or bus. A computer system may havemany additional components such as various adapters or controllers, andserial, parallel and universal bus ports for connection to, e.g.,modems, printers or network interfaces.

In a symmetric multi-processor (SMP) computer, all of the processingunits are generally identical, that is, they all use a common set orsubset of instructions and protocols to operate, and generally have thesame architecture. A typical architecture includes a processor corehaving a plurality of registers and execution units, which carry outprogram instructions in order to operate the computer. The processingunit can also have one or more caches, such as an instruction cache anda data cache, which are implemented using high speed memory devices.Caches are commonly used to temporarily store values that might berepeatedly accessed by a processor, in order to speed up performance byavoiding the longer step of loading the values from a main memorydevice. These caches are referred to as “on-board” when they areintegrally packaged with the processor core on a single integrated chip.

A processing unit can include additional caches, such as a level 2 (L2)cache which may support on-board (level 1) instruction and data caches.An L2 cache acts as an intermediary between the main (system) memory andthe on-board caches, and can store a much larger amount of informationthan the on-board caches, but at a longer access penalty.

A cache has many blocks which individually store the various instructionand data values. The blocks in any cache are divided into groups ofblocks called sets or congruence classes. A set is the collection ofcache blocks that a given memory block can reside in. For any givenmemory block, there is a unique set in the cache that the block can bemapped into, according to preset mapping functions. The number of blocksin a set is referred to as the associativity of the cache, e.g. 2-wayset associative means that for any given memory block there are twoblocks in the cache that the memory block can be mapped into; however,several different blocks in main memory can be mapped to any given set.A 1-way set associative cache is direct mapped, that is, there is onlyone cache block that can contain a particular memory block. A cache issaid to be fully associative if a memory block can occupy any cacheblock, i.e., there is one congruence class, and the address tag is thefull address of the memory block.

An exemplary cache line (block) includes an address tag field, a statebit field, an inclusivity bit field, and a value field for storing theactual instruction or data. The state bit field and inclusivity bitfields are used to maintain cache coherency in a multiprocessor computersystem (to indicate the validity of the value stored in the cache). Theaddress tag is usually a subset of the full address of the correspondingmemory block. A compare match of an incoming address with one of thetags within the address tag field indicates a cache “hit.” Thecollection of all of the address tags in a cache (and sometimes thestate bit and inclusivity bit fields) is referred to as a directory, andthe collection of all of the value fields is the cache entry array.

When all of the blocks in a congruence class for a given cache are fulland that cache receives a request, whether a “read” or “write,” to amemory location that maps into the full congruence class, the cache mustmake one of the blocks in that class available for the new operation.The cache chooses a block by one of a number of means known to thoseskilled in the art (least recently used (LRU), random, pseudo-LRU,etc.). If the data in the chosen block has been modified, that data iswritten to the next lowest level in the memory hierarchy which may beanother cache (in the case of the L1 or on-board cache) or main memory(in the case of an L2 cache). By the principle of inclusion, the lowerlevel of the hierarchy will already have a block available to hold thewritten modified data. If the data in the chosen block has not beenmodified, the value in that block is simply abandoned and not written tothe next lowest level in the hierarchy. This process of freeing up ablock from one level of the cache hierarchy is known as an eviction. Atthe end of this process, the cache no longer holds a copy of the evictedblock. When a device such as the CPU or system bus needs to know if aparticular cache line is located in a given cache, it can perform a“snoop” request to see if the address is in the directory for thatcache.

As microprocessor computing power grows, it becomes more critical forcaches to correspondingly grow in size in order to avoid processingbottlenecks that arise from memory latencies. However, large cachestructures can introduce or exacerbate other problems, such as bandwidthand connectivity. Some high-performance computer systems address theseissues by dividing the cache array and directory into two or moreslices, and allowing multiple access/command ports. One example of sucha sliced cache structure is shown in FIG. 1, which depicts a processingunit 10 having a processor core 12 with on-board instruction and datacaches, and an L2 cache entry array which is divided into two slices 14a and 14 b (slice A and slice B). The L2 cache controller is dividedinto two corresponding slices 16 a, 16 b each having its own directory18 a, 18 b. When processor core 12 issues a load request, the addresstag for the request is sent to one of the directory slices 18 a, 18 b,based on a hash scheme that uses an address bit to direct the request toa given slice (e.g., addr(56)=0 means slice A). The L2 directory sliceperforms the address comparisons and upon detecting a load hit activatesa select signal that controls the output of cache array slices 14 a and14 b. The “addr.rw.ws” signal includes information regarding thecongruence class for the requested memory block, whether the operationis a read or write, and the write set.

Each cache array slice 14 a, 14 b is further divided into four sectors,that is, a given cache line is distributed across all four sectors of aslice. In this example, each cache line is 128 bytes longs, and thedigit pairs in each sector represent the beginning byte number (inhexadecimal) for an 8-byte word of the line, e.g., “00” refers to thefirst 8-byte word in the cache line (bytes 00, 01, 02, 03, 04, 05, 06and 07), and “08” refers to the second 8-byte word in the cache line(bytes 08, 09, 0A, 0B, 0C, 0D, 0E and 0F). Thus, each sector contains 32noncontiguous bytes of a given cache line. All of the sectors are in asingle latency domain but only 32 bytes are output in a given cycle, soit takes four cycles to output a complete 128-byte cache line, with theentire cache array (all sectors) powered up during each of the fourcycles.

Each L2 controller slice 16 a, 16 b has its own read claim (RC), castout (CO) and snoop (SN) machines. Each controller slice further has itsown directory arbiter 20 a, 20 b which handles conflicts between thesemachines and load requests from the processor core. The directoryarbiters are connected respectively to cache arbiters 22 a, 22 b whichcontrol the flow of merge data coming from elsewhere in the memoryhierarchy (e.g., system memory) using separate command ports. Merge flowlogic in each cache slice receives 32 bytes in a given cycle from four8-byte fabric busses that are connected to system memory and variousperipheral devices.

While the use of sliced cache arrays can improve cache bandwidth, thereare still serious problems with power consumption, wiring topology,differential latencies, and recoverability, especially when the designscales to larger cache sizes. As designs grow the cache size by placinglarger numbers of cache array macros, the latency to the farthest arraybecomes multiple clock cycles away from the core compared to the closestcache array. Thus, the prior art mechanism wherein all arrays' accesstimes are in the same clock cycle becomes temporally wasteful, becausethe close arrays must be slowed to match the farthest arrays' accesstime. Although transmission speed can be increased by providing specialwiring (wider/faster), such wiring increases the expense of the designand uses valuable wiring resources, and these problems are compounded indesigns requiring large busses for two cache slices. Even in the case ofa load hit, there can still be a significant delay in accessing andtransmitting the requested cache line, due to the physical layout of thecache and processor core. It would, therefore, be desirable to devise animproved cache structure which could reduce latencies associated with asizeable growth of the cache, particularly latencies arising from loadhits. It would be further advantageous if the cache structure couldmaintain superior directory bandwidth, and still afford a high degree ofrecoverability in the case of a defect in the array.

SUMMARY OF THE INVENTION

It is therefore one object of the present invention to provide animproved cache memory for a computer system.

It is another object of the present invention to provide such a cachememory which is highly scalable to allow large cache arrays withoutsignificantly increasing cache latency.

It is yet another object of the present invention to provide a cachearray topology for large cache structures which takes into considerationdifferent latencies associated with different cache sectors.

The foregoing objects are achieved in a method of operating a cachememory by logically partitioning a cache array having a singleaccess/command port into at least two slices, using a first cachedirectory to access the first cache array slice while using a secondcache directory to access the second cache array slice, and managingaccesses from the first and second cache directories using a singlecache arbiter which controls the single access/command port of the cachearray and manages movement of data inside the cache quadrant and over asingle data bus back to the core. In the illustrative embodiment, thefirst cache directory has a first directory arbiter to handleconflicting access requests within the first cache array slice, thesecond cache directory has a second directory arbiter to handleconflicting access requests within the second cache array slice, and thefirst and second directory arbiters communicate with the cache arbiter.An address tag associated with a load request is transmitted from theprocessor core to either the first or second directory with a designatedbit that associates the address tag with only one of the cache arrayslices whose corresponding directory determines whether the address tagmatches a currently valid cache entry. The cache array may be arrangedwith rows and columns of cache sectors, wherein a given cache line isspread across sectors in different rows and columns, with at least oneportion of the given cache line being located in a first column having afirst latency and another portion of the given cache line being locatedin a second column having a second latency greater than the firstlatency. The cache array outputs different sectors of the given cacheline in successive clock cycles based on the latency of a given sector.

The above as well as additional objectives, features, and advantages ofthe present invention will become apparent in the following detailedwritten description.

BRIEF DESCRIPTION OF THE DRAWINGS

The present invention may be better understood, and its numerousobjects, features, and advantages made apparent to those skilled in theart by referencing the accompanying drawings.

FIG. 1 is a block diagram of a conventional processing unit for acomputer system, depicting a second level (L2) cache memory having asliced directory and array structure;

FIG. 2 is a block diagram of one embodiment of a processing unitconstructed in accordance with the present invention, which includes twoprocessing cores and four L2 cache slice pairs, wherein each processingcore has exclusive access to four of the L2 cache slices;

FIG. 3 is a plan view of a preferred physical layout of the componentsof the processing unit of FIG. 2 as assembled on an integrated circuitchip;

FIG. 4 is a block diagram of an L2 cache memory having a sliceddirectory and unified cache array, constructed in accordance with oneembodiment of the present invention concerned with power conservation;

FIG. 5 is a timing diagram for the data output from the cache array tothe requesting processor core for the cache structure of FIG. 4;

FIG. 6 is a timing diagram for receiving data at the merge flow circuitsof FIG. 4 from the interleaved fabric bus;

FIG. 7 is a schematic diagram of a circuit used for the merge flow ofdata in the L2 cache of FIG. 4; and

FIGS. 8A and 8B together form a block diagrams of another L2 cachememory having a sliced directory and unified cache array, constructed inaccordance with an alternative embodiment of the present inventionconcerned with defect recoverability; and

FIGS. 9A and 9B together form a block diagram of the L2 cache memory ofFIGS. 8A and 8B wherein two interior rows of the cache array have beendisabled due to a defective cache sector.

The use of the same reference symbols in different drawings indicatessimilar or identical items.

DESCRIPTION OF THE PREFERRED EMBODIMENT(S)

With reference now to the figures, and in particular with reference toFIG. 2, there is depicted one embodiment 30 of a processing unitconstructed in accordance with the present invention. Processing unit 30is generally comprised of two processor cores 32 a, 32 b, a sliced level2 (L2) cache 34, two non-cacheable control units (NCUs) 36 a and 36 b(one per core), a fabric bus controller (FBC) 38, and two level 3 (L3)caches 40 a, 40 b with associated L3 controllers 42 a, 42 b. Eachprocessing core 32 a, 32 b includes its own store-through L1 cache(separate program instruction and operand data caches). More than twoprocessing cores may be provided for a single processing unit.

Processing unit 30 may be part of a larger computer system whichincludes various conventional elements (not shown), such as firmware orread-only memory (ROM) and main or random-access memory (RAM) coupled toa peripheral component interconnect (PCI) local bus using a PCI hostbridge. The PCI host bridge can provide a low latency path through whichprocessors 32 a and 32 b may access PCI devices mapped anywhere withinbus memory or I/O address spaces. The PCI host bridge also provides ahigh bandwidth path to allow the PCI devices to access the main memory.Such PCI devices might include a local area network (LAN) adapter, asmall computer system interface (SCSI) adapter providing access to apermanent storage device (e.g., a hard disk drive which stores anoperating system and program files), an expansion bus bridge with userinput devices such as a keyboard and graphical pointer (mouse), an audioadapter, or a graphics adapter. Service processors (not shown) can beconnected to processor cores 32 a, 32 b via a JTAG interface or otherexternal service port, and a processor bridge (not shown) can optionallybe used to interconnect additional processor groups.

The L2 cache portion 34 of the processing subsystem is divided into fourdirectory slice pairs 34 a, 34 b, 34 c and 34 d (for a total of 8slices: A0, B0, C0, D0, A1, B1, C1, D1). Each slice pair has fourmegabytes (MB) of memory, and each processor has exclusive use of fourof the L2 cache slices, i.e., the L2 cache is 8 MB of private memory percore. In the depicted embodiment, core 32 a uses slice pairs 34 a and 34b, and core 32 b uses slice pairs 34 c and 34 d. The individual slicesare selected by an address hashing algorithm residing in cores 32 a, 32b. The same hashing algorithm is used by FBC 38 to route snoop trafficto the L2 slices. For example, bits 55 and 56 of the address can bedecoded to route to slice A when addr(55:56)=00, to slice B whenaddr(55:56)=01, to slice C when addr(55:56)=10, and to slice D whenaddr(55:56)=11.

Each private 4 MB cache is logically partitioned to have two directoryslices and two cache array slices (e.g., A and B). The 4 MB data cacheis further partitioned into four domains or sectors (sect0,1,2,3) wherethe first 32 bytes of a cache line go in sect0 and the last 32 bytes ofthe cache line go in sect3. These sectors are oriented such that allfour sectors sect0,1,2,3 (collectively referred to as quadrant 0) arephysically located together in a corner of the chip, as explainedfurther below. In this example, the cache line size is 128 bytes, L2cache 34 is fully inclusive of the L1 data and instruction caches, andis 8-way set associative. The cache array data and directory array areprotected by error correction code (ECC) having single-bit correctionand double-bit detection (SBC/DBD) capability. A least-recently used(LRU) algorithm is provided and may be enhanced as explained furtherbelow.

L2 slices 34 a, 34 b, 34 c and 34 d generally handle all cacheableload/store instructions, data and instruction prefetches, zero-outinstructions (e.g., the “DCBZ” PowerPC™ instruction), andsynchronization instructions. NCUs 36 a and 36 b handle all othercommunication with the processor cores, such as cache-inhibitedload/store instructions and cache operations (excluding zero-out andsimilar instructions).

FIG. 3 illustrates a general physical layout of the components ofprocessing unit 30 as assembled on an integrated circuit (IC) chip 50 orother substrate. FBC 38 extends horizontally along the central portionof IC chip 50, with core 32 a at the central upper portion and core 32 bopposite the FBC at the central lower portion. Two L2 controllers 44 a,44 b are located between core 32 a and FBC 38 (to service core 32 a),and two more L2 controllers 44 c, 44 d are located between core 32 b andFBC 38 (to service core 32 b). Although the L3 cache array is notincluded as part of the processing unit assembly, the L3 controllers maybe included. In this implementation, L3 controller 42 a-1 interconnectsL2 controller 44 a with L3 cache 40 a; L3 controller 42 a-2interconnects L2 controller 44 b with L3 cache 40 a; L3 controller 42b-1 interconnects L2 controller 44 c with L3 cache 40 b; and L3controller 42 b-2 interconnects L2 controller 44 d with L3 cache 40 b.

The L2 cache array is spread out over four separate quadrants 52 a, 52b, 52 c and 52 d. Each of these quadrants contains the sectors asexplained above, e.g., L2 quadrant 52 a (quad0) contains all foursectors sect0,1,2,3 of a given cache line in L2 slice A0, and alsocontains all four sectors sect0,1,2,3 of a given cache line in L2 sliceB0.

Referring now to FIG. 4, one embodiment of L2 quadrant 52 a is shown ingreater detail (L2 quadrants 52 b, 52 c and 52 d have the samecorresponding structure for the other cache slices C0D0, A1B1, andC1D1). L2 quadrant 52 a has four horizontal cache ways or rows, witheach row containing four sectors (each sector is 16 contiguous bytesarranged in two 4×2-byte arrays). A given row therefore contains 64bytes, or half of a cache line. In the embodiment of FIG. 4, the top andbottom rows together make up cache slice A, and the middle two rowstogether make up cache slice B. The starting sector for a cache line isthe uppermost right array pair together with the lower right array pair(e.g., for cache slice A, the arrays at the far right of the top row andthe far right of the bottom row). The sectors for a cache line progressin order from right to left.

The sectors of L2 quadrant 52 a are arranged in this manner tofacilitate pipelining of the cache output to core 32 a when the cache isfound to contain a currently valid value requested by the core (a loadhit). A load request sent from core 32 a is received by L2 controller 44a, which has a sliced directory. The load address is delivered to eitherthe left (A) or right (B) directory slice based on the setting of adesignated bit in the address field (e.g., addr(56)), and is alsodelivered to a latch that feeds the address to all rows of L2 quadrant52 a. If directory slice A finds a valid matching address, it sends anenable signal (LateSel) to the slice A cache (the top and bottom rows),and if directory slice B finds a valid matching address, it sends anenable signal to the slice B cache (the middle two rows). Each L2directory slice has its own read claim (RC), cast out (CO) and snoop(SN) machines. While the sliced directory effectively has two separateports, the cache is unified with a single access/command port and asingle cache arbiter. The directory arbiters handle access for the RC,CO and SN machines within their respective slices, while the unifiedcache arbiter handles requests for cache access by the RC, CO and SNmachines across both slices. The directory arbiter and cache arbiteralso coordinate their access for certain core requests where thedirectory and cache need access together.

As the address and control information propagates across the cachedomains from right to left, cache controller 44 a selectively enablesthe appropriate cache domains for the read access in successive clockcycles based on their latency. As the cache arrays are read, the datapropagates back from right to left over the reload data bus back to thecore 32 a (multiplexers are used to select between the A and B slicesfor the output). The horizontal (left/right) boundary 54 b betweensectors 1 and 2 thus represents a cycle boundary in this embodiment. Thehorizontal (left/right) boundary 54 a between sector 0 and the corerepresents another cycle boundary.

The cache arbiter sends only one control signal per cycle, but it canschedule different tasks for different sectors to be carried outsimultaneously or in an overlapping manner (such as different stores tobe written to the cache). Only the sectors that are involved with astore operation need to be powered up during a given command sequence.

The timing of the output of L2 cache quadrant 52 a is illustrated inFIG. 5. The core issues the request in cycle 0 with the directoryaddress. Sector0 and sector1 both receive their enable signals andaddress in cycle 1, while sector2 and sector3 see their address in cycle2. The data from the L2 quadrant is delivered to core 32 a on the reloaddata bus at 32 bytes per cycle. The data from all four sectorspropagates back across the L2 quadrant from left to right (crossingthrough the latch boundaries 54 a and 54 b) during cycles 4-6. Duringcycle 5, the data from the arrays labeled 00, 08, 10 and 18 aredelivered to the core. During cycle 6 the data from the arrays labeled20, 28, 30 and 38 are delivered to the core. During cycle 7 the datafrom the arrays labeled 40, 48, 50 and 58 are delivered to the core.During cycle 8 the data from the arrays labeled 60, 68, 70 and 78 aredelivered to the core.

This physical layout allows for pipelining of the cache output in amanner that takes advantage of the differential latencies of the cacheline sectors. Furthermore, by providing a unified cache with outputpipelining, processing unit 30 significantly reduces overall load hitlatency without the need for more expensive cache constructions. Theperformance gain may be further enhanced by using the faster wiresavailable in physical design for the control signal from the core to theL2 controller and then out to the cache arrays as well as for outputfrom the cache to the core. By designating these paths to use wiresconstructed of a premium metal which are up to 16 times the width and 16times the speed of the smallest wires used on the chip, the designerdedicates the fastest wire resource to one of the most sensitiveperformance areas (L2 latency). These features result in a highlyscalable design which keeps load hit latency low in spite of a largegrowth in cache size.

The unified cache structure for two directory slices provides a poweradvantage to the circuits by only having to build the large supportingquad dataflow structure once (as opposed to the prior art structurewhere the dataflow/cache was built once per directory slice). Also, bypartitioning the cache quad into individual sectors, the L2 control forstore operations only needs to enable the cache sector that the store istargeting, as opposed to the prior art which would cause all caches tolight up.

Returning to FIG. 4, if the load or store request from core 32 a missesthe cache (i.e., the cache does not currently contain a valid copy ofthe memory block), then L2 controller 44 a forwards the request tosystem memory via FBC 38. The data is retrieved from elsewhere in thememory hierarchy (e.g., from the L3 cache or system memory) and issubsequently input to the cache in a pipelined fashion, using merge flowcircuits 58 located at the base of cache quad 52 a. Merge flow circuits58 are controlled by the cache arbiter of L2 controller 44 a via a“Mrgflow_ctl” signal. Commands from the cache arbiter to the cache array(to output data) can overlap with commands for merge flow operations.For store requests that are L2 hits, the cache arbiter only activatesthe sector (or sectors) that contains store data. Consequently, twodifferent stores that are directed to different sectors can be performedat the same time.

The read data from FBC 38 is pipelined for a given sector using eight8-byte fabric busses and eight multiplexers. Each multiplexer has twofabric bus inputs, one from interleave bus A (ILVA) and one frominterleave bus B (ILVB). The output of a given multiplexer is connectedto an input line of one of the merge flow circuits 58, e.g., the tworightmost merge flow circuits 58 receive the output of the multiplexerswhose inputs are designated for the first sector of a cache line, i.e.,the first 32 bytes to be stored in the arrays labeled 00, 08, 10 and 18(“ILVA_Byte00_data” and “ILVB_Byte00_data”), while the two leftmostmerge flow circuits 58 receive the output of the multiplexers whoseinputs are designated for the last sector of the cache line, i.e., thelast 32 bytes to be stored in the arrays labeled 60, 68, 70 and 78(“ILVA_Byte60_data” and “ILVB_Byte60_data”). Bus interleaving isscheduled by FBC 38 to avoid conflicts at each of the 8-bytemultiplexers (the notation of “A” and “B” interleave busses is unrelatedto the notation of “A” and “B” for the cache slices).

The timing of the read data for a load or store miss is illustrated inFIG. 6. The address tags for four sectors (e.g., sectors 0, 1, 2, 3)being sent on interleave bus A (“IVLA tag”) are transmitted over a tagbus during the first four cycles of the read operation (i.e., duringcycles 0-3), and the address tags for four sectors being sent oninterleave bus B (“IVLB tag”) are also transmitted over four cycles onanother tag bus, but one cycle behind the interleave bus A tags (i.e.,during cycles 1-4). This tag information is used by L2 cache controller44 a to identify when data is arriving into the cache quad 52 a and bythe merge flow logic 58 to know which data to take. Data transmissionbegins in the third cycle on the IVLA_byte00 data bus, whichsequentially transmits 32 bytes of data (d00 a, d08 a, d10 a and d18 a)over the time span of cycles 2-5. Transmission of the other interleave Adata busses is staggered by one cycle, i.e., the IVLA_byte20 data bussequentially transmits 32 bytes of data (d20 a, d28 a, d30 a and d38 a)over cycles 3-6, the IVLA_byte40 data bus transmits 32 bytes of data(d40 a, d48 a, d50 a and d58 a) over the time span of cycles 4-7, andthe IVLA_byte60 data bus transmits 32 bytes of data (d60 a, d68 a, d70 aand d78 a) over the time span of cycles 5-8. Transmission of theinterleave B data busses is one cycle behind the correspondinginterleave A data bus, i.e., the IVLB_byte00 data bus transmits 32 bytesof data (d00 b, d08 b, d10 b and d18 b) over the time span of cycles3-6, the IVLB_byte20 data bus transmits 32 bytes of data (d20 b, d28 b,d30 b and d38 b) over the time span of cycles 4-7, the IVLB_byte40 databus transmits 32 bytes of data (d40 b, d48 b, d50 b and d58 b) over thetime span of cycles 5-8, and the IVLB_byte60 data bus transmits 32 bytesof data (d60 b, d68 b, d70 b and d78 b) over the time span of cycles6-9.

FIG. 6 shows the relationship of how 32-byte sectors might be scheduledon IVLA and IVLB. The data associated with two different tags (e.g., t00a versus t20 a) may or may not carry data associated with the same128-byte cache line.

Referring now to FIG. 7, the retrieved data from FBC 38 for a load orstore miss is input to a read claim queue (RCQ) array 62 of merge flowcircuit 58. The output of RCQ array 62 is connected to a firstmultiplexer 64 to process store misses. The select line of multiplexer64 is connected to control logic 66 which parses the control signal fromthe L2 cache arbiter to pass the store miss data from RCQ array 62 on toan error-correction code (ECC) generator 68. The output of ECC generator68 is one input to a second multiplexer 70, also controlled by controllogic 66. RCQ array 62 is also connected directly to multiplexer 70 formoving load miss data into the cache array; the ECC generate can bebypassed since the fabric bus has already performed the errorcorrection. Multiplexer 70 then transmits the read data to the cachearrays.

For a store operation, the store data from the core is held in a storequeue array 60, whose output is connected to another input ofmultiplexer 64 and similarly propagates to the cache arrays viamultiplexer 70. Store commands from the cache arbiter to the merge flowcircuits within a given sector can again overlap, i.e., sending a secondstore command while earlier store data is being merged. In addition, thecache arbiter can initiate commands to have different sectors processingdifferent store operations at the same time. For store hits, the cachearbiter only needs to activate the sector(s) that contain the store data(e.g., an 8-byte store to address 00 only needs to affect sector0).Stores can be collected and batch processed by controller 44 a. Thevertical (top/bottom) boundary 56 between the top two rows and thebottom two rows thus represents a cycle boundary for latching the mergeflow data. If the store operation misses the cache and the congruenceclass is already full, a cache line must be evicted. In this case, theevicted cache line passes through an ECC check 72 and acast-out/push/intervention (CPI) array 74 on its way to system memory.

Those skilled in the art will appreciate that the vertical bus structureof the L2 cache quads is generally provided for the merge flowmaintenance (load misses, and stores), while the horizontal busstructure of the L2 cache quads is generally provided for reload (loadhits).

While the processing unit embodiment of FIG. 4 is favored for powerconservation, the present invention further contemplates an embodimentwhich is favored for defect recoverability. This embodiment is shown inFIGS. 8A and 8B, which illustrate one L2 quadrant 52 a′ with itscorresponding controller 44 a′. L2 quadrant 52 a′ still has the samesector arrangement, that is, sector0 being located along the right side(lowest latency), and sector3 being located along the left side (highestlatency), similar to the L2 quadrant 52 a of FIG. 4. However, in theembodiment of FIG. 5, a given row of L2 quadrant 52 a′ is used forportions of both the A and B slices, that is, the set of eight cacheways in a given array are divided across both slices. The top rowcontains selected sectors of the first four members in the congruenceclass for both slices, i.e., A(0:3) and B(0:3), and the bottom rowcontains the remaining sectors of the first four members of both slice Aand slice B. The second row from the top contains selected sectors ofthe last four members in the congruence class for both slices, i.e.,A(4:7) and B(4:7), and the second row from the bottom contains theremaining sectors of the last four members of both slice A and slice B.The command lines from the directory slices in controller 44 a′(LateSel) are accordingly routed to each row of L2 quad 52 a′, e.g., theselect line for the first four members from either directory slice arerouted to the top and bottom rows, while the select line for the lastfour members from either directory slice are routed to the two middlerows.

This construction facilitates partial utilization of the cache (a“half-good” mode) whenever a portion of the cache becomes unusable. Inthe example shown in FIGS. 9A and 9B, at least one of the sectorscorresponding to blocks A(4:7) and B(4:7) has become defective. Theblock(s) may have an original manufacturing defect or a defect thatarose later from, e.g., stray radiation or electrostatic discharge. Thedefect may be discovered by array testing at power-on, or run-timelogging of repeated ECC problems in a given physical array. Once theflaw is discovered, software can flush the L2 cache contents to memory,then direct controller 44 a′ to disable both rows associated with thedefective block. While the example of FIGS. 9A and 9B illustratesdisabling the two middle rows, if the defect had occurred in one of thesectors corresponding to blocks A(0:3) and B(0:3) then the controllerwould disable the top and bottom rows.

In the example of FIG. 9A and 9B two rows are entirely disabled, whichis useful for the situation wherein an address decode circuit is broken,affecting more than one member. For those situations where only a singlemember is defective, a slightly more complicated embodiment could bedesigned to disable only that member and salvage the cache with ⅞capacity.

In this recoverability embodiment, it is preferable to utilize the cacheeviction algorithm to effectuate the disabling of the defective rows.The eviction algorithm may be a least-recently used (LRU) or pseudo-LRUalgorithm which allocates a cache block of the congruence class foreviction based on how recently or frequently that block has beenaccessed. The LRU logic, which resides in controller 44 a′, can includeconfiguration bits or flags set by software to identify defective waysand prevent allocation of any blocks in those ways.

In this manner, if a defect occurs in the cache, the cache can continueoperations with full directory access and all control queues availablewhile only reducing the size of the cache by half (or less), instead ofdisabling an entire cache slice. This approach is particularlybeneficial for certain technical applications which take advantage ofqueue size.

Although the invention has been described with reference to specificembodiments, this description is not meant to be construed in a limitingsense. Various modifications of the disclosed embodiments, as well asalternative embodiments of the invention, will become apparent topersons skilled in the art upon reference to the description of theinvention. For example, while the invention is particularly useful forlarger caches, it is more generally applicable to any type of memoryhierarchy having one or more caches, including non-uniform memory access(NUMA) structures. It is therefore contemplated that such modificationscan be made without departing from the spirit or scope of the presentinvention as defined in the appended claims.

1. A method of operating a cache memory in a given level of a multilevelcache hierarchy, comprising: logically partitioning a cache array of thecache memory into at least first and second slices wherein the firstslice contains a first plurality of bytes arranged in at least a firstrow of the cache array and the second slice contains a second pluralityof bytes arranged in at least a second row of the cache array, the firstplurality of bytes and the second plurality of bytes further beingarranged in columns defining common sectors; receiving a load request ata controller of the cache memory, the load request including an addressfor a requested memory block, and the controller having a firstdirectory associated with the first slice and a second directoryassociated with the second slice; feeding the address to each of therows in the cache array; delivering the address selectively to only oneof the first and second directories based on a setting of a designatedbit in the address; matching the address to an entry in the selected oneof the first and second directories and responsively sending an enablesignal from the selected one of the first and second directories to acorresponding one of the first and second rows; and successivelypowering the sectors of the cache array to pipeline the requested memoryblock to an output of the cache array using a single cache arbiter ofthe cache controller.
 2. The method of claim 1 wherein: the secondplurality of bytes are arranged in the second row and in a third row;and the first plurality of bytes are arranged in the first row and in afourth row.
 3. The method of claim 1 wherein the cache arbiter schedulesdifferent tasks for different sectors to be carried out simultaneously.4. The method of claim 1 wherein the cache arbiter powers differentsectors in successive clock cycles based on the latency of a givensector.
 5. The method of claim 1 wherein: the first directory includes afirst directory arbiter to handle conflicting access requests within thefirst slice; the second directory includes a second directory arbiter tohandle conflicting access requests within the second slice; and thecache arbiter manages access requests from both the first directoryarbiter and the second directory arbiter.
 6. The method of claim 1wherein the cache array includes one or more merge flow circuits foreach sector which receive data to be stored in a designated sector ofone of the rows of the bytes, and the cache arbiter controls the mergeflow circuits to selectively store the data in the designated sector.